1. Field
The present disclosure relates to computer systems and methods in which data resources are shared among data consumers while preserving data integrity and consistency relative to each consumer. More particularly, the disclosure concerns a mutual exclusion mechanism known as “read-copy update.”
2. Description of the Prior Art
By way of background, read-copy update (also known as “RCU”) is a mutual exclusion technique that permits shared data to be accessed for reading without the use of locks, writes to shared memory, memory barriers, atomic instructions, or other computationally expensive synchronization mechanisms, while still permitting the data to be updated (modify, delete, insert, etc.) concurrently. The technique is well suited to both uniprocessor and multiprocessor computing environments wherein the number of read operations (readers) accessing a shared data set is large in comparison to the number of update operations (updaters), and wherein the overhead cost of employing other mutual exclusion techniques (such as locks) for each read operation would be high. By way of example, a network routing table that is updated at most once every few minutes but searched many thousands of times per second is a case where read-side lock acquisition would be quite burdensome.
The read-copy update technique implements data updates in two phases. In the first (initial update) phase, the actual data update is carried out in a manner that temporarily preserves two views of the data being updated. One view is the old (pre-update) data state that is maintained for the benefit of read operations that may have been referencing the data concurrently with the update. The other view is the new (post-update) data state that is seen by operations that access the data following the update. In the second (deferred update) phase, the old data state is removed following a “grace period” that is long enough to ensure that the first group of read operations will no longer maintain references to the pre-update data. The second-phase update operation typically comprises freeing a stale data element to reclaim its memory. In certain RCU implementations, the second-phase update operation may comprise something else, such as changing an operational state according to the first-phase update.
FIGS. 1A-1D illustrate the use of read-copy update to modify a data element B in a group of data elements A, B and C. The data elements A, B, and C are arranged in a singly-linked list that is traversed in acyclic fashion, with each element containing a pointer to a next element in the list (or a NULL pointer for the last element) in addition to storing some item of data. A global pointer (not shown) is assumed to point to data element A, the first member of the list. Persons skilled in the art will appreciate that the data elements A, B and C can be implemented using any of a variety of conventional programming constructs, including but not limited to, data structures defined by C-language “struct” variables. Moreover, the list itself is a type of data structure.
It is assumed that the data element list of FIGS. 1A-1D is traversed (without locking) by multiple readers and occasionally updated by updaters that delete, insert or modify data elements in the list. In FIG. 1A, the data element B is being referenced by a reader r1, as shown by the vertical arrow below the data element. In FIG. 1B, an updater u1 wishes to update the linked list by modifying data element B. Instead of simply updating this data element without regard to the fact that r1 is referencing it (which might crash r1), u1 preserves B while generating an updated version thereof (shown in FIG. 1C as data element B′) and inserting it into the linked list. This is done by u1 acquiring an appropriate lock (to exclude other updaters), allocating new memory for B′, copying the contents of B to B′, modifying B′ as needed, updating the pointer from A to B so that it points to B′, and releasing the lock. In current versions of the Linux® kernel, pointer updates performed by updaters can be implemented using the rcu_assign_pointer( ) primitive. As an alternative to locking during the update operation, other techniques such as non-blocking synchronization or a designated update thread could be used to serialize data updates. All subsequent (post update) readers that traverse the linked list, such as the reader r2, will see the effect of the update operation by encountering B′ as they dereference B's pointer. On the other hand, the old reader r1 will be unaffected because the original version of B and its pointer to C are retained. Although r1 will now be reading stale data, there are many cases where this can be tolerated, such as when data elements track the state of components external to the computer system (e.g., network connectivity) and must tolerate old data because of communication delays. In current versions of the Linux® kernel, pointer dereferences performed by readers can be implemented using the rcu_dereference( ) primitive.
At some subsequent time following the update, r1 will have continued its traversal of the linked list and moved its reference off of B. In addition, there will be a time at which no other reader process is entitled to access B. It is at this point, representing an expiration of the grace period referred to above, that u1 can free B, as shown in FIG. 1D.
FIGS. 2A-2C illustrate the use of read-copy update to delete a data element B in a singly-linked list of data elements A, B and C. As shown in FIG. 2A, a reader r1 is assumed be currently referencing B and an updater u1 wishes to delete B. As shown in FIG. 2B, the updater u1 updates the pointer from A to B so that A now points to C. In this way, r1 is not disturbed but a subsequent reader r2 sees the effect of the deletion. As shown in FIG. 2C, r1 will subsequently move its reference off of B, allowing B to be freed following the expiration of a grace period.
In the context of the read-copy update mechanism, a grace period represents the point at which all running tasks (e.g., processes, threads or other work) having access to a data element guarded by read-copy update have passed through a “quiescent state” in which they can no longer maintain references to the data element, assert locks thereon, or make any assumptions about data element state. By convention, for operating system kernel code paths, a context switch, an idle loop, and user mode execution all represent quiescent states for any given CPU running non-preemptible code (as can other operations that will not be listed here). The reason for this is that a non-preemptible kernel will always complete a particular operation (e.g., servicing a system call while running in process context) prior to a context switch.
In FIG. 3, four tasks 0, 1, 2, and 3 running on four separate CPUs are shown to pass periodically through quiescent states (represented by the double vertical bars). The grace period (shown by the dotted vertical lines) encompasses the time frame in which all four tasks that began before the start of the grace period have passed through one quiescent state. If the four tasks 0, 1, 2, and 3 were reader tasks traversing the linked lists of FIGS. 1A-1D or FIGS. 2A-2C, none of these tasks having reference to the old data element B prior to the grace period could maintain a reference thereto following the grace period. All post grace period searches conducted by these tasks would bypass B by following the updated pointers created by the updater.
Grace periods may be synchronous or asynchronous. According to the synchronous technique, an updater performs the first phase update operation, then blocks (waits) until a grace period has completed, and a second phase update operation, such as removing stale data, is performed. According to the asynchronous technique, an updater performs the first phase update operation, specifies the second phase update operation as a callback, then resumes other processing with the knowledge that the callback will eventually be processed at the end of a grace period. Advantageously, callbacks requested by one or more updaters can be batched (e.g., on a multi-part callback lists) and processed as a group at the end of an asynchronous grace period. This allows asynchronous grace period overhead to be amortized over plural deferred update operations. An RCU state machine mechanism is used to start and end grace periods and advance the RCU callbacks, one grace period at a time, through several stages of callback list processing. A typical RCU callback list may comprise three or four sublist portions that segregate the batched RCU callbacks into callback groups that are processed at the end of different grace periods. A callback must advance through each sublist before it is ready for invocation. In RCU implementations used in the Linux® kernel, there is typically one RCU callback list per processor, and list tail pointers are used to divide each callback list into four sublists. The callback sublists are named for their corresponding tail pointers. Newly arrived callbacks that must await a subsequent grace period before they can be invoked are placed on a first sublist named “RCU_NEXT_TAIL.” As grace periods elapse, the callbacks first advance to a second sublist named “RCU_NEXT_READY_TAIL,” then to a third sublist named “RCU_WAIT_TAIL,” and finally arrive at a sublist named “RCU_DONE_TAIL.” All callbacks on the RCU_DONE_TAIL sublist are deemed to be ready for invocation.
More recently, RCU grace period processing has been adapted to account for processor low power states (such as, on Intel® processors, the C1E halt state, or the C2 or deeper halt states). Operating systems can take advantage of low power state capabilities by using mechanisms that withhold regular timer interrupts from processors (in a low power state) unless the processors need to wake up to perform work. The dynamic tick framework (also called “dyntick” or “nohz”) in existing versions of the Linux® kernel is one such mechanism. In RCU implementations designed for low power applications in the Linux® kernel, a compiler configuration option called RCU_FAST_NO_HZ is available. This option allows processors to be placed in low power states even if there are pending RCU callbacks, provided none require immediate invocation and the processor is not needed for grace period advancement processing. Such processors will awaken after a short interval (e.g., four scheduling clock periods), at which time the processor will attempt to advance its RCU callbacks. This capability results in significant power savings for some workloads.
Unfortunately, the RCU_FAST_NO_HZ option can also result in greatly increased grace period latencies. This is due to the fact that the processors which are sleeping with callbacks cannot take full advantage of subsequent grace periods. A waking processor reacts only to a change in a tracked grace period completion number. It does not account for the actual number of additional grace periods that have elapsed. So even if several grace periods elapse while the processor was sleeping, the processor will take advantage of only one, thus potentially delaying its callbacks for another sleep period. On a busy system, a callback will normally take roughly 1.5 grace periods to advance through the callback sublists. After arriving on the RCU_NEXT_TAIL sublist, the callback will advance to the RCU_NEXT_READY_TAIL sublist and as soon as the callback becomes known to the RCU subsystem. When the next grace period starts, the callback advances to the RCU_WAIT_TAIL sublist. When that grace period ends, the callback advances to the RCU_DONE_TAIL sublist for invocation. In the case of a sleeping processor, if the processor goes idle before the callback advances to the WAIT sublist, an additional 0.5 grace periods will be required to advance the callback when the processor wakes up. This is because the processor does not receive credit for more than one of the potentially many grace periods that elapsed while it was idle. As a consequence, the scheduling clock tick is typically not deactivated for processors that have callbacks.
Another scenario causing increased grace period latency for a sleeping processor (in a RCU_FAST_NO_HZ kernel) is when no other processor in the system needs a grace period to start. In that case, the start of the next grace period will be delayed until the sleeping processor awakens, further degrading grace period latency for another sleep period. Because no other processor sees any reason to start a new grace period, the RCU subsystem remains idle while the processor is asleep. When the processor wakes up, an additional grace period is required to advance its callbacks. Had the RCU subsystem been aware of the processor's need for an additional grace period while the processor slept, the processor could have instead woken up to find that the grace period that it needed had already completed. Because of this latency issue, the RCU_FAST_NO_HZ option causes processors that are sleeping with callbacks to attempt several times to advance the RCU state machine on the off-chance that some of the callbacks can then be invoked. This state machine work often has no effect and consumes processor time, and thus energy.
Accordingly, there is a need for a technique that allows processors to sleep with RCU callbacks without needing to expend large amounts of processor time on idle entry, while taking full advantage of any grace periods that elapse during the sleep interval, and by initiating grace periods as needed during the sleep interval without having to awaken to do so.